A well-structured framework for analysing petri net extensions

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  • Information and Computation 195 (2004) 129

    www.elsevier.com/locate/ic

    A well-structured framework for analysing petri net extensions

    Alain Finkela ,, Pierre McKenzieb ,1 , Claudine PicaronnyaaLaboratoire Spcication et Vrication (CNRS URA 8643), cole Normale Suprieure de Cachan, 61, Avenue du

    Prsident Wilson, 94235 Cachan Cedex, FrancebInformatique et recherche oprationnelle, Universit de Montral, C.P. 6128, succ. Centre-Ville, Montral,

    Que., Canada H3C 3J7

    Received 4 December 1998; revised 23 May 2003

    Abstract

    Transition systems dened from recursive functions p p are introduced and named WSNs, or well-structured nets. Such nets sit conveniently between Petri net extensions and general transition systems. In therst part of this paper, we study decidability properties of WSN classes obtained by imposing natural restric-tions on their dening functions, with respect to termination, coverability, and four variants of the bounded-ness problem. We are able to precisely answer almost all the questions which arise, thus gaining much insightinto old and new generalized Petri net decidability results. In the second part, we specialize our analysis toWSNsdened fromafne functions, which elegantly encompassmost Petri net extensions studied in the litera-ture.Again,we studydecidability properties of natural classes of afneWSNwith respect to the above six com-putational problems. In particular, we develop an algorithm computing limits of iterated nonnegative afnefunctions, in order to decide the path-place variant of the boundedness problem for non-negative afneWSN. 2004 Published by Elsevier Inc.

    Keywords: Logic and verication of computer program; Theory of parallel and distributed computation; Algorithms;Decidability

    Corresponding author. Fax: +33 1 47 40 75 29.E-mail address: nkel@lsv.ens-cachan.fr (A. Finkel).1 Research performed in part while on leave atUniversitt Tbingen. Supported by the (German)DFG, the (Canadian)

    NSERC, and the (Qubec) FCAR.

    0890-5401/$ - see front matter 2004 Published by Elsevier Inc.doi:10.1016/j.ic.2004.01.005

  • 2 A. Finkel et al. / Information and Computation 195 (2004) 129

    1. Introduction

    A transition system is a set S endowed with a transition relation , i.e., with a binary re-lation on the set S . It is a WSTS, i.e., a well-structured transition system, if a reexive and tran-sitive well-quasi-ordering on S exists which is compatible with (i.e., such that s1 t1and s1 s2 imply the existence of a sequence of transitions leading from t1 to t2 s2 [9,1,11]). Wedeal in this paper with innite state transition systems, with WSTS, and with Petri net extensions[3,4,14,17,20].

    Our motivating theme is that Petri net extensions can naturally and fruitfully be dened andstudied from the perspective of WSTS. This theme is only marginally new: WSTS were introducedas abstractions of Petri nets in the rst place, as a means of identifying the essential properties ofPetri net problems, algorithms, and extensions. But WSTS in fact capture many other innite statesystems [11], and WSTS are so general that (1) only rather coarse Petri net extension decidabilityresults, concerning boundedness detection, for example, can be dealt with from WSTS, and (2)the classication of Petri net extensions afforded by WSTS is not as complete as one might havehoped.

    In this paper, we respond to both weaknesses of theWSTS with the introduction of theWSN, orwell-structured net. AWSN is aWSTS inwhich the set S isp and the transition relation is giv-en by a nite set of recursive functions with upward-closed domains, i.e., with domainsK p suchthat u K and u v imply v K . We establish the suitability of the WSN model for a systematicstudy of Petri net extensions and their decidability properties.

    Our paper is formed of two parts. The rst part deals with WSN in general. The second partdeals with WSN dened from afne functions.

    In the rst part, we dene progressively stronger properties of functions used in the denitions ofWSNs: nondecreasing (i.e., u < v f(u) f(v)), increasing (i.e., u < v f(u) < f(v)), -increas-ing (i.e., increasing, with the added requirement that knowledge of the components along whichu < v provides information on the components along which f(u) < f(v)), strongly increasing (i.e.,increasing componentwise), each with or without the -recursivity hypothesis that the extensionof the function to the domain ( {})p is recursive. These properties give rise to two groups offour WSN classes, and we exhaustively study these classes with respect to six decision problems.These problems are the usual termination problem (is the reachability tree of a given state s0 nite?),coverability problem (is there a state s in the reachability set of a given state s0 which covers a givenstate s?), boundedness problem (is the reachability set of a given state s0 nite?), place-boundednessproblem (is the component i bounded, given a start state s0?), together with two new variants ofthe boundedness problem introduced here to demonstrate the subtleties of boundedness detection.Let us note (s0) the state reached from s0 by applying . These two variants are the path-un-bounded-witness problem (given s0 and a sequence of transitions such that s0 < (s0), is there anunbounded component in the sequence of states starting at s0 and obtained by repeatedly applying?) and the path-place-boundedness problem (given s0 and a sequence of transitions such thats0 < (s0), is the component i unbounded in the sequence of states starting at s0 and obtained byrepeatedly applying ?).

    Our results are summarized on Figs. 1 and 2: only four of the 48 relevant decidability questions(for some of them, answers were already implicitly or explicitly known) remain open. We also claimnew results on -recursive functions.

  • A. Finkel et al. / Information and Computation 195 (2004) 129 3

    Fig. 1. Abbreviations.

    Fig. 2. Decidability (D) or undecidability (U) for well-structured nets.

    We single out, from the rst part of our paper, the following results:

    Coverability for-WSN, i.e.,WSNwhose functions satisfy the-recursivity hypothesis, is showndecidable, while coverability is shownundecidable for ourmost restrictiveWSNclass (i.e., strong-ly increasing WSN) in the absence of the -recursivity hypothesis.

    The decidability status of the four boundedness problems for WSN is resolved. This provablyseparates three of our WSN classes, and the fourth class is distinguished from the others whenviewed as a -WSN class.

    Boundedness is shown decidable for increasing WSN but undecidable for -WSN. Path-unbounded-witness is shown decidable for -increasing WSN but undecidable for increas-

    ing WSN. Path-place-boundedness, known to be decidable for reset Petri nets (and for transfer Petri nets),

    is shown undecidable for strongly increasing WSN and for -WSN. Place-boundedness is showndecidable for strongly increasing-WSN. This is proved by showing

    that the coverability tree algorithm of Karp and Miller [10,15,16,18] still applies when the exactcomputation of the limit of an innite repetitive sequence is replaced by the computation of agood enough approximation. Yet, the place-boundedness is shown undecidable for both stronglyincreasing WSN and -increasing -WSN.

    In the second part of our paper, we specialize our study to afne WSN, i.e., to WSN denedfrom afne functions. We do this because testing for nontrivial properties of general recursive

  • 4 A. Finkel et al. / Information and Computation 195 (2004) 129

    functions is undecidable, and because, in the concrete cases likely to be encountered in practice,functions are either afne, or their properties are known a priori, or these are easy to test. In Fig.3, we record some easy (but appealing) connections between afne WSN, Petri net extensions, andgeneral WSN. For example, generalized transfer Petri nets can be seen as WSN dened from afnefunctions f(X) = AX + B satisfying A 0 with every column of the matrix A different from 0.Afne WSN thus capture a wealth of known Petri net extensions, including Petri nets themselves[23], Reset Petri nets [7], Post SM nets [27], Double Petri nets, and they suggest new extensions ofindependent interest.

    We then developed a nontrivial algorithm to solve the following problem: given a nonnegativep p-matrix A and two nonnegative p-vectors B and X , compute d 1 such that limn f nd (X)exists, where f is the afne function f(X) = AX + B, and compute this limit. We use this algorithmto solve the path-place-boundedness problem for afne WSN dened from nonnegative matricesand vectors. This WSN class is incomparable with Petri Nets and it does not have an immediatepractical value but it provides an interesting instance in which the path-place boundedness problemis decidable.

    Finally, we record the decidability status of our six computational problems, for ve prominentclasses of afne WSN, including Petri nets and Self-Modifying nets [27,28]: only two out of the30 resulting questions (whose answers were known for Petri nets and Self-Modifying nets) remainopen. Fig. 4 reports these results.

    Section 2 in this paper introduces our various functions, the notion of limit, -recursivity, andproves relevant undecidability results. Section 3 formally denes well-structured nets and our com-putational problems. Section 4 examines decidability of termination and coverability for WSN.Section 5 does the same for the four boundedness problems, Section 6 summarizes our decidability

    Fig. 3. The world of well-structured nets. An arrow C D indicates C D and a double line indicates equality.

  • A. Finkel et al. / Information and Computation 195 (2004) 129 5

    Fig. 4. Afne -well-structured nets.

    results for WSN. Section 7 discusses afne WSN, their relationship with WSN and extended Petrinets, and their decidability properties. Section 8 concludes.

    2. Notation and preliminaries

    Sets.When X is a set, a word v X is a possibly empty nite sequence of elements of X ; we write|v| for the length of v.

    We write [p] for {1, 2, . . . , p}, + for \ {0}, and for the standard completion {} of theset of nonnegative integers (where n < for any n ).Upward-closed sets. Fix any integer p 1. Let be the usual ordering on . We also write forthe ordering on p dened by: for all u, v p, u v if for every i [p], u(i) v(i), where u(i) isthe ith component of u. We write u < v if u v and u /= v. Addition in and in p is denedcomponentwise, with x + = + x = for any x . Let U + V denote the sum of two subsetsU , V p: U + V = {u+ v : u U , v V }. We say that a set K p is upward-closed (for ) ifK = K + p .

    A basis of an upward-closed set K is a set B such that K = B + p . We dene the downward-clo-sure of u p, written u, as u = {v p : v u}. The two well-known facts which follow can befound in [12], or in [11] where their proofs are briey recalled:

    1. Every upward-closed set K p has a nite basis.2. Any strictly increasing sequence K0 K1 K2 of upward-closed sets in p is nite.

    We will require another result due to Valk and Jantzen:

    Theorem 2.1 ([30]). A nite basis of an upward-closed set K p is effectively computable iff for anyvector u p the predicate u K /= is decidable.

    Transition systems.A transition system (TS) is a structure S = S ,, . . .where S = {s, t, . . .} is a setof states, and S S is a set of transitions. For every s S , wewrite Succ(s) (respectively Pred(s))for the set {t S : s t} of immediate successors of s (respectively for the set {t S : t s} of im-mediate predecessors of s). We write s

    + t to mean that there exist k > 0 and a sequence of statess0 = s, s1, . . . , sk = t such that s0 s1 sk . We write s t to mean that s + t or s = t. Wesay that S is nitely branching if Succ(s) is nite for each s S; the usefulness of this property istied to Knigs lemma [19] stating that an innite tree, nitely branching, has an innite path.

  • 6 A. Finkel et al. / Information and Computation 195 (2004) 129

    Reachability. A state t of a TS S is reachable from a state s if s t.

    The reachability set of S from s0 is denoted RS(S , s0) and is dened as the set of states reachablefrom s0.

    The reachability tree RT(S , s0) of S with the initial state s0 is a directed unordered tree wherenodes are labelled by the states of S . The root node is labelled by s0. A node labelled by s has ason labelled by t iff t Succ(s). Moreover, any two different sons of a node must have differentlabels.

    Orders and limits. Consider an innite nondecreasing sequence (un)n 1 of natural numbers. Thissequence is bounded if there exists v such that un v holds for every n . We dene

    limn un =

    { if (un)n 1 is not bounded,max{un : n 1} otherwise.

    Now consider a nondecreasing innite sequence (un)n 1 of p-tuples un p . We also dene:

    limn un =

    (limn un(1), limn un(2), . . . , limn un(p)

    ).

    For our purpose, we shall need the following orders: for any nonempty subset Q [p], let

  • A. Finkel et al. / Information and Computation 195 (2004) 129 7

    gf(x) = g(f(x)). The notion of growth signature will help us dene classes of functions which willillustrate the subtleties of the boundedness question:

    Growth signature. The growth signature of an arbitrary function f : p p is the function fdened as follows:

    f : {Q [p] : Q /= } 2{Q[p]:Q /=}Q {Q [p] : u, v Dom f , [u

  • 8 A. Finkel et al. / Information and Computation 195 (2004) 129

    (3) If f and g are recursive, then gf is recursive.(4) If f is of type t, 1 t 4, then f is of type s,1 s t.

    Proof. Straightforward. We nonetheless argue (1) for the case of -increasing functions. Let f andg be -increasing functions. Consider /= Q [p]. Pick u, v Dom gf . Then u, v Dom f andf(u), f(v) Dom g. Hence by the -increasing properties of f and g,

    u

  • A. Finkel et al. / Information and Computation 195 (2004) 129 9

    g(m,) = g(m 1,)+{(2, 0) if the TMm halts,(1, 0) otherwise.

    The rest of the section studies the impact of the properties of f on the decidability properties of f .Consider adapting in the natural way the nondecreasing, increasing, and strongly increasing

    properties to the case of functions f : p for p > 1. Then, increasing and strongly increasingbecome synonymous properties. Furthermore, as for any u in p having at least one componentequal to , we must have f (u) = , the extension f of any increasing total recursive functionf : p is recursive. On the other hand, projecting onto its rst component the function g usedin proving Proposition 2.4 yields a nondecreasing total recursive function g : 2 that is not-recursive.

    Returning to nondecreasing total recursive functions f : , whose extensions f are recur-sive as we have seen, we might hope to be able to compute the description of a Turing machinecomputing f , given one computing a nondecreasing total f (this is clearly possible when f isincreasing). The next proposition shows that this is impossible. Let NDEC be the set of Turing ma-chinesm computing a nondecreasing total recursive function fm : , and let NDEC NDECcomprise those m for which fm() = . Now assume to the contrary, that there is a total recursivefunction g such that g(m) is a machine computing f m whenever m NDEC . Let L be the languageof all m such that the machine g(m) halts on and outputs . Then L is a recursively enumerablelanguage such that L NDEC = NDEC, which contradicts the following:

    Proposition 2.5. Any language L such that

    L NDEC = NDECis not recursively enumerable.

    Proof. Consider such an L. Note that L is not necessarily an index set so that Rices theorem (forrecursively enumerable index sets) does not apply. Consider for each m the Turing machine M(m)computing the function f(m) : dened as

    f(m)(n) ={t if the TMm halts in exactly t steps and t < n,n otherwise.

    Note that(m) NDEC for everym. Now theTMm does not halt iff f (m)() = iff(m) NDECiff (m) L NDEC iff (m) L. Hence, the recursive map is a many-one reduction fromthe complement of the halting problem to L, proving that the latter is not recursivelyenumerable.

    2.2. Transitions as afne functions

    An afne2 function f : p p is given by f(X) = AX + B for some A = (Ai,j) pp andB p and a recursive denition domain Dom f {X p : AX + B 0}.

    2 We restrict the usual sense of afne functions. In our context, all afne functions have a non negative linear part.

  • 10 A. Finkel et al. / Information and Computation 195 (2004) 129

    We allowDom f to differ from {X p : AX + B 0}, to be able to simulate transitions of Petrinets. For instance, suppose that, in a Petri net with p places, the transition t only tests whether thereis a token in place 1. This transition is the identity function t(X) = X (hence A = Id and B = 0) andt(X) 0 for everyX p . Yet its denition domain isDom t = + p1 and it is not equal top .

    We denote by Aj the jth column of the matrix A, for any j in [p]. We write A 0 to mean thatAi,j 0, for all i, j and A Id to mean that for all i,Aii 1 and A 0.

    The following holds for every afne function f dened by f(X) = AX + B:

    1. f is nondecreasing,2. f is increasing iff f is -increasing iff Aj /= 0, for every j in [p],3. f is strongly increasing iff A Id .

    Hence, given an afne function f bywayof its deningmatrixA and vectorB, it is straightforwardto determine its increasing property. Moreover, belonging to the set {X p : AX + B 0} iseasily tested, and it is easy to see that this set is upward-closed whenever f is nondecreasing (A 0).

    The composition of two afne functions is afne and every afne function is recursive. Further-more, the extension of a nondecreasing afne function f is recursive as the extension can also bedened by its afne form AX + B on X p with the usual conventions 0 = 0 and m = , forall m such that m > 0.

    3. Well-structured nets

    We recall [11] that awell-structured transition system (WSTS) is a transition system S = D,,equipped with a reexive and transitive relation D D fullling the following two conditions:

    (1) well-quasi-ordering: for any innite sequence s0, s1, . . . D, there exist two indices i < j withsi sj .

    (2) compatibility: is (upward) compatible with, i.e., for any s1 t1 and transition s1 s2, thereexists t2 such that t1

    t2 and s2 t2.

    Well-structured nets. A well-structured net (WSN) is a triple S = p , F , for some p , where F isa nite set of nondecreasing recursive functions dened on upward-closed subsets of p .

    Intuitively, an -well-structured net is a generalization of a WSN dened on p. -Well-struc-tured nets.A -well-structured net (-WSN) is a triple S = p, F , for some p , where F is a niteset of nondecreasing recursive functions dened on upward-closed subsets ofp such that for everyfunction f F , we have f = fNf where fNf : Nf Nf is the restriction of f on Nf .A WSN S = p , F , or a -WSN S = p, F , are, respectively,

    increasing if every function in F is increasing, -increasing if every function f F is -increasing and an algorithmwhich computes the growth

    signature of f is given, strongly increasing if every function in F is strongly increasing.

  • A. Finkel et al. / Information and Computation 195 (2004) 129 11

    Note that a WSN S = p , F , gives rise to the WSTS p , F,, where s Ft for (s, t) p p is dened to mean that f(s) = t for some f F . Therefore, we use for S all notions de-ned for p , F, (reachability...). In the context of a WSN, a place is an integer; we can speak ofthe boundedness of a place: given S and an initial state s0, the place i is bounded if there exists k such that for each s RS(S , s0), s(i) k . Note also that everyWSN is nitely branching, because F isnite.

    Remark 3.1.Every Petri netN with p place is aWSN S = p , F , such that every function (calledtransition) t F may be written as t(X) = X + vt , where vt is an integer vector of dimension p . ResetPetri nets are extended Petri nets in which a transitionmay clear (reset) a place. Formally, a reset Pe-tri net N with p places is aWSN such that every function t in F is an afne function t(X) = AtX + vtin which the matrix (of integers) At satises Id At 0, where Id and 0 are the identity and thenull matrices respectively, and inequality is taken componentwise.

    Let S = p, F , a-WSN, s1, s2 p and = fkfk1 . . . f1, with f1, ..., fk F . We write s1 s2to mean that there exist t1, ..., tk+1 such that s1 = t1 f1t2 f2t3 fktk+1 = s2.

    Given a state so in p and = fkfk1 . . . f1, with f1, ..., fk F , we dene RS(S , s0, ) = {s p | s0

    n2 s, n , 2 prex of }.

    Proposition 3.2 ([9]). Let S = p , F , be a WSN, s1 and s2 two states in p , and a nite sequencesuch that s1

    s2 and s1 s2. Then there is a sequence of states (sn)n 1 such that for all n, sn sn+1and sn sn+1.

    Problems. We will consider the following computational problems, where in each case, a WSNS = p , F , or a -WSN S = p, F , is given as part of the input:

    1. Termination. Given a state s0 in p , is the reachability tree RT(S , s0) nite?2. Coverability. Given states s0 and s inp , does there exist a state s in RS(S , s0) which covers s, i.e.,

    which satises s s?3. Boundedness. Given a state s0 in p , is the reachability set RS(S , s0) nite?4. Path-unbounded-witness. This problem is not a decision problem, but rather a search problem in

    the sense of [22]: given a state s0 inp and a nite non empty sequence such that s0(s0) and

    s0 < (s0), determine if there is an unbounded place inRS(S , , s0) (i.e., a place i satisfying k ,s RS(S , , s0) such that s(i) k) and if so identify3 at least one such an unbounded place i.

    5. Path-place-boundedness: Given a state s0 inp , a sequence such that s0(s0) and s0 < (s0),

    and a place i, is the place i bounded in RS(S , , s0) (i.e., does there exist k such that s(i) k ,s RS(S , , s0))?

    6. Place-boundedness: Given an initial state s0 and a place i, is i bounded?

    3 Although this problem is not a decision problem, we say that the problem is decidable if there exists a Turingmachinewhich solves it (and always halts).

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    In this list, we have introduced two new problems related to the boundedness problem. ThePath-place-boundedness can be seen as the place-boundedness problem restricted to a particularpath of the reachability tree. It may be decidable when the boundedness problem is not (e.g.reset Petri Nets [7]). Note that decidability of the path-place-boundedness problem implies de-cidability of the path-unbounded-witness problem. As we show later, the reverse implication isfalse.

    We do not mention the usual Reachability problem (given two states s and s, does s shold?), as it is irrelevant to our paper, being undecidable in all the eight Petri Net extensions weshall study (cf. Fig. 1).

    4. When are termination and coverability decidable?

    For every state s we denote by s the set {t D : t s}.

    Theorem 4.1. Termination is decidable for well-structured nets.

    Proof. Let S = p , F , be a WSN. If s1 fs2 for a function f F and s1 t1, then becausef is nondecreasing and Dom f = Dom f , we have t1 ft2 and s2 t2. This property and therecursivity of f allow us to apply Theorem 4.6 from [11] and then to decide termination.

    Theorem 4.2. Coverability is decidable 4 for -well-structured nets.

    Proof.We will show that it is possible to compute a nite basis of Pred(s) for any s p . Thenby using Theorem 3.6 from [11], we will conclude that coverability is decidable. Fix s p and writePredf (s) for {t p | f(t) s}. We have:

    Pred(s) =fF

    Predf (s) =fF

    Predf (s) =fF

    Predf (s),

    where the rightmost equality uses the fact that Predf (s) = Predf (s), which holds becauseDom f = Dom f and f is nondecreasing, for every f F . Because F is nite, it sufces to beable to compute a nite basis of Predf (s) for f F . By Theorem 2.1, namely the Valk and Jantzenresult, a nite basis of Predf (s) is computable iff the predicate t Predf (s) /= is computablefor all t p. And this holds because

    t Predf (s) /= iff (t p , t t)[f(t) s] iff f (t) s,where the left to right implication in the second iffrelies on the non decreasing property of f .Now the inequality f (t) s can be checked because f is recursive.

    4 Such a statement means: there is a TM deciding the coverability problem, given the premise that the input WSN isan -WSN.

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    The computability of f is a hypothesis needed to decide coverability, since we have as a coun-terpart to Theorem 4.2 that our strongly increasing WSNs have an undecidable coverability whenthis hypothesis is absent:

    Theorem 4.3. The coverability problem for strongly increasing well-structured nets is undecidable.5

    Proof. Consider the family {fj : 2 2} of strongly increasing recursive functions

    fj(n, k) ={(n, 0) if k = 0 and TMj runs for more than n steps(n, n+ k) otherwise

    and the strongly increasing function g : 2 2 dened by g(n, k) = (n+ 1, k). The strongly in-creasing WSN Sj = 2, {fj , g}, with initial state (0, 0) has the property that the state (1, 1) iscoverable iff the TMj halts. Hence there is no Turing machine which correctly determines cover-ability and halts whenever its input is a strongly increasing WSN.

    5. When are boundedness problems decidable?

    5.1. The boundedness problem

    An increasing WSN veries the property that if s1 s2 and t1 > s1, then there exists t2 such thatt1 t2 with t2 > s2, while we can only conclude t2 s2 in the case of a WSN. Boundedness forincreasing WSNs can thus be decided by searching for a sequence s0

    s1 s2 with s1 < s2. Then, can be iterated to produce an increasing sequence of reachable states s0

    s1 s2 s3 . . . withs1 < s2 < s3 . . ., and the system is unbounded. Because anyWSN by denition has a recursive Succ(i.e., Succ(s) is computable, for every s in p ), Theorem 4.11 from [11] yields:

    Theorem 5.1. Boundedness is decidable for increasing well-structured nets.

    We note further that Theorem 5.1 is tight. If we omit the increasing property from the hypoth-esis of Theorem 5.1, and yet add the -recursivity property, the boundedness problem becomesundecidable:

    Theorem 5.2. Boundedness is undecidable for -well-structured nets.

    Proof. Boundedness is undecidable for reset Petri nets [7], which are afne WSNs and every afneWSN is a -WSN (see Section 2.2).

    5 The precise technical meaning of this statement is that there is no Turing machine which correctly decides whethersome state in RS(S , s0) covers s, and halts, whenever its input S is indeed a strongly increasing WSN (with its deningfunctions given by Turing machines) together with s0 and s.

  • 14 A. Finkel et al. / Information and Computation 195 (2004) 129

    5.2. The path-unbounded-witness problem

    In this section, we show how to detect at least one unbounded place in some unbounded WSNs.Indeed we show the problem decidable for -increasing WSNs. This decidability result provablydoes not extend to increasing WSNs, as we also prove in this section.

    Theorem 5.3. The path-unbounded-witness problem is decidable for -increasing WSN.

    Proof. Consider a -increasing WSN S = p , F ,. Consider the edge-labelled graphG = ({Q [p] : Q /= }, {(Q, f ,Q) : Q f (Q)}),

    where f is the (computable) growth signature of f . Consider any non empty nite path in thisgraph from a node Q to a node Q. Let f1, f2, . . . , fk be the sequence of edge labels encounteredalong the path. A straigthforward induction on k proves that

    u, v Dom fkfk1 f1, [u

  • A. Finkel et al. / Information and Computation 195 (2004) 129 15

    This implies that all the places in the non-empty set Q grow unbounded when is iterated from s0.Hence any place in Q is a satisfactory answer.

    Theorem 5.3 is tight. It provably extends neither to the guaranteed identication of at least oneunbounded place in increasing well-structured nets, as we now show, nor to the decidability ofpath-place-boundedness in -increasing WSN, as we will show in the next section.

    Theorem 5.4. The path-unbounded-witness problem is undecidable for increasing well-structured nets.

    Proof. For each m, we dene f m : 2 2 by

    f m(k , n) ={(k , n+ 1) if the TMm halts in at most k 1 steps(k + 1, n) otherwise.

    Every f m is total (hence upward-closed), recursive, and increasing.Now, from any initial state (k , n) 2, the increasingWSN Sm = 2, {f m}, possesses one and

    only one unbounded place, which is the second place iff the TMm halts. Suppose to the contrary thatthere is a TM M which solves the path-unbounded-witness problem. On input s1 = (0, 0), = f m,and s2 = (1, 0),M will necessarily come up with the unique unbounded place resulting from the iter-ation of f m from s1. This place is the second place if TMm halts, and the rst place otherwise, whichmeans thatM solves the halting problem. Hence no suchM exists, and so path-unbounded-witnessproblem is undecidable.

    5.3. The path-place-boundedness problem

    Theorem 5.5. Path-place-boundedness is undecidable for strongly increasing well-structured nets andfor -well-structured nets.

    Proof.

    Consider the strongly increasing WSN Sj = 2, {fj , g}, from the proof of Theorem 4.3, withinitial state (0, 0) and the sequence = gfj . Then, place 2 is unbounded along the innite path(0, 0)

    ... if and only if TMj halts. Consider the family of functions gj : dened by

    gj(n) ={n if TMj halts in at most n stepsn+ 1 otherwise.

    For every j 1, the function gj is nondecreasing, total, and -recursive (gj() = ). LetTj = , {gj},with initial state 0. The only place is bounded along the innite path 0 gj ... gj ...if and only if the TMj halts.

    The following proposition justies our interest in the path-place boundedness problem.Althoughboundedness in undecidable for Reset Petri nets [7], path-place boundedness is decidable.

  • 16 A. Finkel et al. / Information and Computation 195 (2004) 129

    Proposition 5.6. Path-place-boundedness is decidable for reset Petri nets.

    Proof . Given a reset Petri net N with p places, a marking m, a place i and a sequence of transi-tions such that m

    m1 with m < m1. By monotonicity, ( is non-decreasing), the sequence maybe innitely repeated; let us write the innite path m

    m1 m2 .Let us write = fkfk1...f1. Because each afne function fi(X) = AiX + vi satises 0 Ai Id

    and vi is a vector of integers, we deduce that there exist A and v such that is the afne function(X) = AX + v with still 0 A Id and v a vector of integers (where A is computable fromthe Ais and v from the vis and Ais).

    For every n and 1 j p , we have: if does not contain any reset operation on j (i.e.,A(j, j) = 1)then mn(j) = m(j)+ nv(j); moreover, if contains a reset operation on place j, then the sequence(mn(j))n 1 is stationary and for every n, 2 n, mn(j) = m1(j) and hence place j is bounded alongthe path m

    m1 m2 .We will now prove that i is unbounded along the innite path m

    m1 m2 iff m1 m2and m1(i) < m2(i).

    (1) If i is unbounded along the innite path m m1 m2 , then cannot contain any reset

    operation on place i (hence A(i, i) = 1), v must be positive (otherwise the path would be nite);hence, we obtain that for every n, 1 n, mn mn+1. Moreover, we can nd two integers r andk such that: 1 r, 1 k , mr mr+k , and mr(i) < mr+k(i).Hence, frommr(i) < mr+k(i) and 0 < k , we obtain 0 < v(i). Hence, we also have:m1(i) < m2(i).

    (2) Let us prove that if m1 m2 and m1(i) < m2(i) then i is unbounded along the innite pathm

    m1 m2 . First, by monotonicity, ( is non-decreasing), the path can be innite-ly continued and then 0 v . Because m1(i) < m2(i), the sequence cannot contain any re-set operation on place i and then 1 v(i). From mn(i) = m(i)+ nv(i), we obtain that mn(i)goes to innity when n goes to innity. Hence, the place i is not bounded on the innite pathm

    m1 m2 .

    5.4. The place-boundedness problem

    Let us recall the construction of the coverability tree of a Petri net, which is used to solve theplace-boundedness problem. The strategy of the construction is to develop every path of the reach-ability tree until one meets two markings m and m such that m0

    m m with m

  • A. Finkel et al. / Information and Computation 195 (2004) 129 17

    To apply the above strategy to construct a coverability tree for a WSN, it is rst necessary tobe able to effectively extend functions to p. This means that functions have to be -recursive.Furthermore, the extension f of any nondecreasing function f is nondecreasing, but any type-iproperty of f may not be preserved on f . In particular, strongly increasing functions extend-ed to p may no longer be increasing (for example, when f is dened as f(m, n) = (m,m+ n),f(, 0) = f(, 1)).

    Moreover, either limn n(u)must be computable (but this is not true for nondecreasing -re-cursive functions), or a good enough approximation lof limn n(u)must be computable such thatu < l limn n(u) and l has at least one more -component than u, for every innite increasingsequence (n(u))n 1.

    Let -max be the following function which takes two elements s1 and s2 in p and returns an

    element in p, dened by:

    function -max(s1, s2): returns s : ps s2for each i such thats1(i) < s2(i) and s2(i) /= dos(i) ;

    We will only apply this function to pairs of states (s1, s2) satisfying s1 s2. Note that in gen-eral, -max can satisfy -max(u, (u)) < limn n(u), for example when f(m, n) = (m+ n, n+1) and u = (0, 0), in which case f(u) = (0, 1) and -max(u, f(u)) = (0,), while limn f n(u) =(,).

    Nevertheless, any strongly increasing WSN satises the interesting property:

    Proposition 5.7. Let p , F , be a strongly increasing WSN and let u, v be in p. If there existf1, . . . , fk in F such that u

    v for = f k f k1 . . . f 1, and there exists a nonempty Q such thatu v, and -max(u, v) contains |Q| more s than u.

    Proof.As (n(u))n 1 is a nondecreasing sequence, limn n(u) exists. As a composition of stronglyincreasing functions, is strongly increasing. So u

    v and u

  • 18 A. Finkel et al. / Information and Computation 195 (2004) 129

    ly (n1, s1)+KMT (n, s)) for saying that there is path labelled by (respectively a possibly empty,

    respectively non-empty) path in KMT from node (n1, s1) to the node (n, s).

    KarpMiller_tree (p, F , : -WSN, s0 : state in p) returns KMT: tree

    NODES ; ARCS ;UNPROCESSED_NODES {(0, s0)}; {* Initial singleton set of unprocessed nodes,each node referred to as (node number, temporary label) *}while a node (n, s) is in UNPROCESSED_NODES do

    UNPROCESSED_NODES UNPROCESSED_NODES \ {(n, s)};for all nodes (ni, si) such that (ni, si)

    KMT (n, s) and si < s dos -max(si, s);

    NODES NODES {(n, s)}; {* Add node and fix label forever *}if no node (n1, s) exists such that (n1, s)

    +KMT (n, s) thenfor each function f F such that s Dom f do

    n new node number;UNPROCESSED_NODES UNPROCESSED_NODES {(n, f (s))};ARCS ARCS {(n, f , n)};

    KMT (NODES ,ARCS);Before proving that the algorithm terminates on all -WSNs, we remark some useful elementary

    facts about the nite paths of the KarpMiller tree.

    Lemma 5.8.LetS=p, F ,bean-WSN, s0 an initial state inp andKMT(S , s0) theKarpMillertree of (S , s0).We have:

    If (0, s0) KMT (n1, s) KMT (n2, s) then the node n2 has no successor(along a branch of the KMT, the same state cannot appear more than twice).

    if (n, s) KMT (n, s) then (s) s. if (n, s) KMT (n, s) and s < s then s has at least one more than s.

    Proposition 5.9. The KarpMiller tree algorithm terminates when it is applied to an -WSN.

    Proof. Suppose that the KarpMiller tree algorithm does not terminate and so KMT is innite.Because the set of functions is nite, the innite tree KMT is nitely branching and there is aninnite branch in KMT:

    (0, s0) KMT (n1, s1) KMT .... KMT (nk , sk) KMT ...Because is a well-ordering on p, we may extract an innite non-decreasing sequence {ski} thatcannot be stationary (by Lemma 5.8).

    So we may also choose an innite increasing sequence along this branch:

    (nk0 , sk0)KMT (nk1 , sk1) KMT .... with ski < ski+1

  • A. Finkel et al. / Information and Computation 195 (2004) 129 19

    We recall that (n, s)KMT (n, s) and s < s imply that smust contain at least onemore than

    s, by Lemma 5.8.After a nite number (at least p) of such executions, no new can be inserted. Hence, for every

    n, skp = skp+n and the sequence (skn)n 1 is stationary, which produces a contradiction. Hence KMTis nite.

    In the following, we will say a state s in KMT(S , s0), for a state s such that there exists a node(n, s) in KMT(S , s0).

    When the KarpMiller algorithm is applied to an -WSN, it produces a nite tree such thatevery reachable state is smaller than (or equal to) a state in KMT (see Proposition 5.11). But ingeneral, the KMTmay very roughly cover the reachability set, in the following sense: some state inthe KMT may be strictly greater than the limit of any innite sequence of reachable states. Whenthe KarpMiller algorithm is applied to a strongly increasing -WSN, it produces a tree, oftencalled a coverability tree, which nely covers the reachability set: each state in KMT is smallerthan (or equal to) the limit of an innite sequence of reachable states. The following property ofstrongly increasing functions (with upward-closed domain) is used for proving that the cover isne.

    Lemma 5.10. Let f : p p be any strongly increasing function with upward-closed domain. Let(un)n 1 be a nondecreasing sequence in Dom f and u its limit in

    p. There exists a strictly increasing

    sequence of natural numbers ((n))n 1 such that limn f n(u) = limn f n(u(n)).

    Proof. Let us recall that any nondecreasing sequence in either goes to , either isstationary. For all m, we dene Im as the set of places i such that the nondecreasing sequence(f m(un)(i))n 1 is non stationary. Let i Im. Then for all integer n, there exists p > n such thatf m(un)(i) < f

    m(up)(i). Because f is strongly increasing, f m+1(un)(i) < f m+1(up )(i). Therefore i Im+1. We have shown the sequence (Im)m1 is nondecreasing (for ), therefore there exist a sub-set I of places and a natural number N such that Im = I for all m N . These I and Nsatisfy:

    (1) For i I andm N , the sequence (f m(un)(i))n 1 is non stationary, thus its limit value is . Wemay write:m N , 7m > 0 such that i I , n 7m, we have f m(un)(i) > m.

    (2) For i / I and m N , the sequence (f m(un)(i))n 1 is stationary. Thus the limit value must befm(u)(i). In that case, we may write:

    m N , 8m > 0 such that i / I , n 8m, we have f m(un)(i) = f m(u)(i).

    For all m N , we dene by induction (m) as the maximum of Am, Bm, and (m 1)+ 1. Wethen have:

    (1) For i I and m N , f m(u(m))(i) > m.(2) For i / I and m N , f m(u(m))(i) = f m(u)(i).

    Hence, we have limn fn(u) = limn f n(u(n)).

  • 20 A. Finkel et al. / Information and Computation 195 (2004) 129

    Proposition 5.11.For every -WSN S = p, F , with initial state s0, we have:

    1. for every state s RS(S , s0), there exists a state s KMT(S , s0) such that s s,2. If, moreover S is strongly increasing, we have: for every state s KMT(S , s0), there exists an

    innite non decreasing sequence (sn)n 1 of reachable states sn RS(S , s0) such that s limn sn.

    Proof .

    1. Let s be a reachable state in RS(S , s0); we will use induction on the length of with s0

    s. When is the empty word, the statement holds because it is always true that s0 KMT(S , s0).Consider a sequence (in the reachability tree) f of length k + 1. We dene inductively f = f .From s0

    s f f(s) and the induction hypothesis, we deduce that there exists (0, s0) KMT(n, s) with s p and s s. Because s s and f is nondecreasing, we obtain f (s) f (s).Now there are two possible cases:

    The node (n, s) has no successor. This is necessarily because there exists a path (n1, s) KMT(n, s).Because f (s) exists, the node (n1, s) has at least one successor (n, t) such that: (n1, s)

    fKMT(n, t) with f (s) t.

    The node (n, s) has at least one successor then there is an arc labelled f : (n, s) fKMT (n, t)with f (s) t.

    In both cases, we have: there exists a state t in KMT such that t f (s) f(s).2. We use induction on the number of times k one has used the function -max on a sub-path (in

    KMT) from the initial state to state s.

    Let k = 0; suppose that there is a branch (0, s0) KMT (n, s) such that the function -maxhas not been used along this branch. Then s is a reachable state and we may write sn = s forevery n.

    Now let k = n+ 1, and consider (0, s0) KMT (n1, s) KMT (n2, t) in which the lastapplication (in the KarpMiller algorithm) of the function -max occurred directlybefore n2, whence: t = -max(s, (s)). Then by the induction hypothesis, there exists aninnite nondecreasing sequence of reachable states (sn)n 1 such that s limnsn;and from the algorithm, we may deduce that s

  • A. Finkel et al. / Information and Computation 195 (2004) 129 21

    Remark 5.12. From the last proposition, we may prove that for every innite nondecreasing se-quenceof reachable states (sn)n 1, there exists a state t in theKarpMiller tree such that limnsn t;in general, there is no reason for having: limnsn = t. Moreover, every state t in KMT isnot necessarily the limit of an innite nondecreasing sequence of reachable states. Hence forstrongly increasing -WSN, the KarpMiller tree does not exactly enjoy the same properties asthose dened for Petri nets. But we still have a strong relation between the KarpMiller treeand the reachability set: KMT(S , s0) =RS(S , s0) where KMT(S , s0) is the downward-closureof the set of states occurring in KMT(S , s0); this is sufcient for deciding the place-boundednessproblem.

    We thus obtain:

    Theorem 5.13. For a strongly increasing -well-structured net p, F ,, place-boundedness isdecidable.

    Proof.FromProposition 5.11, we show that a place i is not bounded in S = p, F ,with the initialstate s0 iff there is anode (n, t) inKMT(S , s0) such that t(i) = . Bydenition, i is notbounded in (S , s0)means that there exists an innite sequence of reachable states (tn)n 1 such that limntn(i) = .Hence, from Proposition 5.11 (1), we obtain that if i is not bounded in (S , s0) then there existst KMT(S , s0) such that t(i) = . Conversely: suppose there is a t KMT(S , s0) such that t(i) = .FromProposition 5.11 (2), there exists an innite non-decreasing sequence of reachable states (sn)n 1such that: t limnsn. Hence, limnsn(i) = and i is not bounded in (S , s0). Because the KMTalgorithm terminates (cf. Proposition 5.9) for-well-structured nets, we conclude that place-bound-edness is decidable.

    Theorem 5.13 is optimal in the following sense:

    Theorem 5.14. Place boundedness is undecidable for both strongly increasing well-structured nets andfor -increasing -well-structured nets.

    Proof. Undecidability of place-boundedness holds for -increasing -WSN because transferPetri nets [7] are -increasing -WSN and undecidability of place-boundedness holds forthose. For proving the undecidability for strongly increasing WSN, we again use the family{Sj} with initial state (0, 0) as in Theorem 4.3. The place 2 is unbounded if and only if TMjhalts.

    6. Summary

    Recall the six problems dened in Section 3. We have shown in Sections 4 and 5 the followingresults:

    Theorem 6.1.The decidability status of the six problems dened in Section 3 of this paper, for the eightclasses depicted in Fig. 1, is summarized on Fig. 3.

  • 22 A. Finkel et al. / Information and Computation 195 (2004) 129

    7. Afne realizations of well-structured nets

    7.1. On the difculty of testing general recursive functions

    Given a set F of functions p p prescribed by their respective Turing machines, decidingwhether F gives rise to a WSN is undecidable, a consequence of Rices theorem [25] which we stateas Proposition 7.1:

    Proposition 7.1. For each type (t) of function chosen from (1) nondecreasing, (2) increasing, (3)-in-creasing, and (4) strongly increasing, the following languages

    (a) {k | the kth TM computes f : of type (t)}(b) {k | the kth TM computes f : of type (t) whose domain is upward-closed}(c) {k | the kth TM computes f : whose domain is upward-closed}

    are undecidable.

    Nevertheless, in concrete cases, it is generally known or easy to decide in which class the giv-en set of functions lies. In many Petri net extensions, this given set is in fact contained in theset of afne functions. Many connections between afne functions and Petri net extensions areshown in Fig. 2. Section 7.2 describes WSN for which deciding the properties of relevant func-tions is easy. Section 7.4 summarizes the decidability status of problems involving afnefunctions.

    7.2. Petri net extensions and afne well-structured nets

    In Section 2.2, we discussed the fact that every afne WSN is a -WSN.Figs. 3 and 4 locate Petri net extensions within our classes of afne functions. Fig. 3 also locates

    the afne classes within our general framework of well-structured nets.The Double Petri nets [7] appearing on Fig. 3 are extended Petri nets in which some arc from a

    transition to a place p may double the content of p . Double Petri nets are then afneWSN such thatevery afne function f(X) = AX + B satises Id A 2Id . Transfer Petri nets [7] are extended Petrinets in which some transitions may transfer the content of a place p to another place p . TransferPetri nets are then afne WSN such that every afne function f(X) = AX + B satises A 0 andevery column Aj satises :iAij = 1. Generalized Transfer Petri nets may transfer and duplicate, sothat they are afne WSN such that A 0 and every column Aj is nonnull. Reset Petri nets [2,7] areextended Petri nets in which an arc may clear a place. This class corresponds to afne WSN suchthat every afne function f(X) = AX + B satises Id A 0.

    Many Petri net extensions are thus surprisingly well parametrized by afne functions.

    7.3. An algorithm for path-place-boundedness

    A nonnegative afne function is an afne function f : p p such that f(X) = AX + B withB 0.

  • A. Finkel et al. / Information and Computation 195 (2004) 129 23

    A nonnegative afne WSN is a WSN such that every function f F is nonnegative afne.6 Weprove here that path-place-boundedness is decidable for nonnegative afne WSN.

    For f a nonnegative afne function and X a nonnegative vector inp , we prove that the limit off nd (X) for a suitable d is computable.We can determine preciselywhich entries in f nd (X) go to inn-ity and thereforewhich entries in f nd+r(X).7 go to innity, for all r such that 0 r d 1. Further-more, we show that this limit is computable by searching a computable graph. Let f(X) = AX + Bbe a nonnegative afne function. Note that, for n , f n is also a non negative afne function,precisely: f n(X) = AnX + (An1 + An2 + + A+ I)B.

    Recall that multiplication is extended by 0 = 0 and x = , x {0}. Lets rst remarkthe following:

    Lemma 7.2. Let (An)n 0 be a sequence of non negative matrices having a limit A with possibly coefcients. Let X be any non negative vector. Then limn AnX = AX.

    So we may restrict the problem to the study of the sequences (An)n 0 and (I + A+ + An)n 0.

    Notations. Let A be a matrix. We write A = 0 or A = when each coefcient of A is respectively 0or . We write A for limn An when the limit exists.

    Let A = (Ai,j)(i,j)[p][p] be a matrix with nonnegative integers coefcients. Let us denote An =(A

    (n)i,j ) (where the brackets in the exponent are used to avoid confusion with the (i, j)-entry of A

    raised to the nth power).We associate to A a directed graphGA whose vertices are 1, . . . , p andwhose set of edges is dened

    by: there is an edge leaving vertex i and entering vertex j if and only if Ai,j /= 0; then, the coefcientAi,j is called the weight of this edge. For a subgraph C of GA, we shall denote by AC the submatrix(Ai,j)(i,j)CC . We shall need the general decomposition of a non-negative matrix A in term of itsassociate graph GA, as it is usually done in the study of nite homogeneous Markov chains or theproof of Perron-Frobenius theorem (cf. [13], [21], [26]).

    A path of length n from a vertex i to a vertex j is a sequence i = i0, i1, . . . , in1, in = j such that(ik ,ik+1) is an edge inGA, for k=0, 1, . . . , n 1.Theweightof this path is theproductAi,i1Ai1,i2 . . . Ain1,j .The path i = i0, i1, . . . , in1, in = j is called a cycle whenever n /= 0 and i = j. The cycle i = i0, i1, . . . ,in1, in = i is called a circuit if i0, i1, . . . , in1 are distinct. The index of imprimitivity of a graph is theg.c.d. of the lengths of all circuits of the graph. When this index of imprimitivity is equal to 1, wesay the graph is primitive. We rst recall two properties of this index of imprimitivity:

    Lemma 7.3. Let G be a graph whose index of imprimitivity is denoted by d. Let i be a vertex of G.Then all cycles going through i in G have lengths multiple of d.

    6 A nonnegative afne function f : X AX + B in a WSN satises A 0 and B 0, as f is required to be nonde-creasing.

    7 If we denote by L(X) the limit of (f nd (X))n 0, whenever d > 1, the d sub-sequences (f nd (X))n 0,(f nd+1(X))n 0, . . . , (f nd+d1(X))n 0 are convergent with respective limits L(X), f(L(X)), . . . , f d1(L(X)).

  • 24 A. Finkel et al. / Information and Computation 195 (2004) 129

    Proof. Let i = i0, i1, . . . , in1, in = i be a cycle of length n going through i. Let us show that n is amultiple of d by induction: If i0, i1, . . . , in1 are distinct, then this cycle is a circuit and n is a multipleof d by denition. Otherwise, let k , l with 0 k < l n be such that ik = il. Take l as small aspossible. So ik , ik+1, . . . , il is a circuit of length l k , which implies l k is a multiple of d . Andi0, i1, ik , il+1 . . . , in1 is a cycle of length n l+ k < n, thus a multiple of d by induction.

    Lemma 7.4. Let G be a strongly connected graph whose index of imprimitivity is denoted by d. Let ibe a vertex of G. Then d is the g.c.d. of the lengths of all cycles going through i in G.

    Proof. Let e be the g.c.d. of the lengths of all cycles going through i in G. Thanks to Lemma 7.3e is a multiple of d . Let j = j0, j1, . . . , jn1, jn = j be a circuit in G. Let i = i0, i1, . . . , il = j be apath going from i to j and let j = il+1, il+2, . . . , im = i be a path going from j to i in G. Then, i =i0, i1, . . . , il = j0, j1, . . . , jn = il+1, il+2, . . . , im = i and i = i0, i1, . . . , il = il+1, il+2, . . . , im = i are cyclesgoing through i, thus their lengths are multiple of e. thus, the length of the circuit is a multiple of e.This shows d is a multiple of e.

    As we shall see, the asymptotic of the sequence (An)n 0 may be completely described by thegraph GA. The proof is based on the fact that A

    (n)i,j is the sum of the weights of all paths of length n

    from i to j.The next two results correspond to particular cases of the nal result of this section.

    Lemma 7.5. Let C be a strongly connected component of GA whose index of imprimitivity is equalto 1. Then, either C is reduced to a singleton and AC is one of the 1 1 matrices (0) or (1), orAC = .

    Proof . Recall some facts about primitive strongly connected graphs : We discard the case wherethe graph C has no edge at all, i.e., C is a singleton and AC = (0). Let i, j in C . Note that becauseC is strongly connected there exists a path from i to j (so there exists t such that A(t)i,j > 0) andthere exists a path from i to j which goes through all vertices of C . Because C is maximal for thisproperty, all paths from i to j in GA are in fact in C . Furthermore, we may nd circuits in C withlengths l1,l2,. . ., lk such that g.c.d.(l1, l2, . . . , lk) = 1. With a path of length l0 from i to j crossing allthese circuits, we can build a path in C of length l0 + n1l1 + n2l2 + + nklk , for any non negativeintegers n1, n2, . . . nk . It means there exists a path from i to j of any large enough length, i.e., q such that A(n)i,j > 0 for n > q.

    If A(n)i,i = 1, for sufciently large integers n, then there is only one circuit in AC with all edgeslabelled by 1. So AC is a permutation matrix, which means GAC is a disjoint union of circuits. ButGAC = C is strongly connected and the index of imprimitivity of C is equal to 1, so AC = (1).

    Otherwise, there exists s such that A(s)i,i 2. Let i, i1, . . . , ir1, j a path of length r from i to jand let q be such that A(n)j,j > 0, if n q. Now, if n s+ q+ r, we may decompose n as u+r + ts with u {q, . . . , q+ s 1} such that u n r modulo s and t > nqsrs . Then A(n)i,j (A

    (s)i,i )

    tAi,i1 . . . Air1,jA(u)j,j 2t which shows limn A

    (n)i,j = .

  • A. Finkel et al. / Information and Computation 195 (2004) 129 25

    Lemma 7.6.LetC be a strongly connected component ofGA with index of imprimitivity equal to c > 1.Let D,E two strongly connected components of GAcC and (i, j) D E. Then

    (Ac)i,j =0 if D /= E1 if {i} = D = E = {j} and A(c)i,i = 1 otherwise.

    Proof. Let us rst remark that there exists a path of length l from i to j in GAcC if and only if thereexists a path of length cl from i to j in GA. Let D be a strongly connected component of GAcC and ibe a vertex of D. Thanks to Lemma 7.4, the index of imprimitivity of D is the g.c.d. of the lengths ofall cycles in GAcC going through i, i.e., the quotient by c of the g.c.d. of the lengths of all cycles in GAgoing through i. Therefore, it is equal to 1. LetD,E two strongly connected components ofGAcC and(i, j) D E. If D /= E, there cannot exist a path from i to j and a path from j to i in GAcC . Let ussuppose, for example, that i cannot be joined to j by a path in GAcC , which means i cannot be joinedto j by a path of length multiple of c in GA. As all cycles have lengths multiple of c, j cannot bejoined either to i by a path of length multiple of c. So n, A(nc)i,j = A(nc)j,i = 0. When D = E, we applyLemma 7.5 to (A(c)(i,j))(i,j)DD which denes a strongly connected graph with index of imprimitivityequal to 1.

    The general case: The next theorem explains how to describe the behaviour of the sequence(An)n 0, without making any assumption of strong connectivity or on the index of imprimitivityof the graph GA associated to the matrix A. We compute the strongly connected components ofGA. Grouping nodes by components, we nd a permutation matrix P such that PAP1 is a triangu-lar-blocks matrix where each diagonal block is irreducible (corresponding to a strongly connectedcomponent of GA). Let d be the l.c.m. of the indices of imprimitivity of all components of GA. Wealready know by the two previous lemma how to compute the limit of the powers of each irreduciblediagonal block of Ad . Let r be the number of strongly connected components of GAd . We partitionthe matrix B = PAdP1 = (Bi,j)(i,j)[r][r] such that each block Bi,i corresponds to a strongly con-nected component of GAd . We shall call this component a 0, 1 or block according to the nature ofits limit in Lemma 7.5.

    We consider the graph GB: vertices are the strongly connected components of GAd , and edgesare dened by: there is an edge leaving vertex i and entering vertex j if and only if Bi,j /= 0. As thematrix B is upper-triangular, each path i, i1, , it , j satises i i1 it j. We say that sucha path is increasing. We say this path is strictly increasing if i < i1 < < it < j. If 1 i < j r,we denote by Pi,j the set of strictly increasing paths from i to j; this set contains only paths oflength j i so is nite and computable. We may decompose Pi,j = jit=1P ti,j where P ti,j is the set ofpaths of length t in Pi,j , for all t such that 1 t j i. All these sets allow us to express the powersof B:

    B(n)i,j =

    min(ji,n)t=1

    (i,i1,...,it1,j)P ti,j

    p0++pt=nt

    Bp0i,i Bi,i1B

    p1i1,i1 . . . Bit1,jB

    ptj,j.

  • 26 A. Finkel et al. / Information and Computation 195 (2004) 129

    So we get the next result which describe the behaviour of the sequence (And )n 0:

    Theorem 7.7. Let i < j r.

    If there exists a strictly increasing path in GB connecting i to j which goes through a node in a block or through two nodes in different 1 blocks, then Bi,j = . If all strictly increasing paths in GB connecting i to j go through nodes in 0 blocks or if there existsno path connecting i to j, then Bi,j = 0. Otherwise, all strictly increasing paths in GB connecting i to j go through nodes in 0 blocks exceptone 1 block Bl,l. Then Bi,j =

    (i,i1,...,l,...,it1,j)Pi,j Bi,i1Bi1,i2 . . . B

    l,l . . . Bit1,j.

    The sameconsiderations allow to computeL= limn(And + A(n1)d + + I). If limn A(nd)i,j /=0, the i, j-term in L is . If limn A(nd)i,j = 0, the number of paths from i to j is nite ; in particular,they do not cross any cycle and theymust be of length p . So the i, j-term in L is 1+ Adi,j + + Akdi,jwhere k is such that kd p (k + 1)d if i /= j or 1 if i = j.

    Remark 7.8. This result may in fact be extended to matrices with non negative real coefcients(see [24]).

    Theorem 7.9. Let f(X) = AX + B be a nonnegative afne function fromp top . Let X0 be a nonneg-ative vector. Then there exists an integer d 1 such that limn f nd (X0) exists in p. Furthermore,d and limn f nd (X0) are computable.

    Proof . Let us write the informal algorithm which computes d and the limit limn f nd (X0) :

    1. Get the strongly connected components (C) of GA and their imprimitivity indices (dC)2. Compute CdC

    for each irreducible diagonal block of A, corresponding to a strongly connected

    component of GA(i) Search for connected components (E) of CdC(ii) If one of E = (1), then E = (1) for all E(iii) Else, E = for all E

    3. Compute the index d and Ad(using Theorem 7.7)

    4. Compute

    n And

    5. Compute (Ad)X0 + (n And )B (using Lemma 7.2)

    Remark 7.10. Even if its interest seems to be more theoretical than practical, we may very roughlystudy the complexity of this algorithm, in the dimension p of the matrix A and in the size of thecoefcients of A, B, X0, denoted respectively by m(A), m(B), m(X0). Computing strongly connect-ed components [5] and indices of imprimitivity of strongly connected components in GA may bedone by inspecting the matrix A and its ith powers (for i p), thus is polynomial in p . Step 3 inthe algorithm may be most expensive as we have to compute the dth-power of A, where d is thel.c.m. of the indices of imprimitivity of strongly connected components. As all these indices canbe bounded by p , we roughly bound d by p ! The cost of such a computation is therefore bound-

  • A. Finkel et al. / Information and Computation 195 (2004) 129 27

    ed by p !O(p3) in operations and p !O(p3)m(A) in the size of coefcients. The last steps are lessexpensive.

    Corollary 7.11. Path-place boundedness is decidable for nonnegative afne WSN.

    Proof.Using notations of theorem 7.9, the sequence (f n(X0))n 0 is bounded on the ith-componentif and only if all vectors limn f nd (X0), limn f nd+1(X0),. . ., limn f nd+d1(X0) have a niteith-component.

    Interestingly, although the procedure in Theorem 7.9 computes limits and always terminates,we do not see how to bound the number of limit computations required by the coverability treestrategy from Section 5.

    Note that this method cannot be applied if we do not suppose B 0 because of Lemma 7.2, inwhich the hypothesis X 0 is crucial. We could give partial results with other kinds of hypotheses(AB = B which captures the Petri net model, {An | n } is a nite set which captures the transferPetri net model. . .).

    7.4. Afne well-structured nets

    Theorem 7.12. The decidability status of the six problems dened in Section 3 of this paper, for veclasses of afne -WSN, is summarized on Fig. 4.

    Proof.First column: termination and coverability are decidable because afneWSNwith A 0 are-WSN (i.e., in S1); then see Fig. 2. Path-place boundedness, and hence path-unbounded-witness,are decidable because path-place boundedness for afneWSNwith A 0 reduces to the path-placeboundedness for afneWSN with A 0 and B 0: let us write (X) = AX + B and ; = (s0) s0;we have ; = As0 + B s0 0. Now, n(s0) = s0 + ; (1+ A+ A2 + + An1). Let us note fthe afne function f(X) = AX + ; We remark that for every n, f n (0) = ; (1+ A+ A2 + +An1) = n(s0) s0; hence place i is bounded along the innite path obtained in iterating fromthe initial state s0 iff i is bounded along the innite path obtained in iterating f from the initialstate 0. We now remark that the afne function f satises A 0 and ; 0; hence the path-placeboundedness is decidable for f , so we may decide whether i is bounded on f n (0); hence we mayalso decide the path-place boundedness on n(s0) too.

    Boundedness, and hence place-boundedness, are undecidable because boundedness is alreadyundecidable for reset Petri nets [7] and reset Petri nets are afne WSN with 0 A Id ; then seeFig. 3.

    Second column: path-place-boundedness, hence path-unbounded-witness, is decidable (Corol-lary 7.11). The case of boundedness is open; boundedness does not reduce to the path-place-bound-edness problem for (A 0,B 0)-afneWSN directly because one could imagine detecting a nitepath s

    s1 s2 with s1 < s2 and limn sn p .Third column: boundedness is decidable because afne WSN with each column of each matrix

    greater than 0 are increasing -WSN (i.e., in S2).Fourth column: follows from [2729].Fifth column: afne -WSN with A = Id are Petri nets!

  • 28 A. Finkel et al. / Information and Computation 195 (2004) 129

    8. Conclusion

    Wehave answered all but six decidability questions concerning termination, coverability, bound-edness, path-unbounded-witness, path-place boundedness, and place-boundedness, for our eightclasses of general WSN, and for ve classes of afne WSN. These results are summarized on Figs.2 and 4. We highlight specic contributions:

    The coverability problem (and the termination problem) are shown decidable for every -WSN.This uses a characterization of effective upward-closed subsets of p .

    To clarify the delicate role of model variations on the difculty of deciding boundedness, weintroduced two new boundedness problems, called the path-unbounded-witness and the path-place-boundedness problems, as companions to the boundedness and the place-boundednessproblems classically studied. Our WSN classes provably distinguish the decidability propertiesof these four boundedness problem variants.

    The place-boundedness problem remains decidable for any class of WSN in which each functionis -recursive and increases in the strong sense.

    Reset Petri nets, Generalized transfer Petri nets, transfer Petri nets, post-SM nets, double Petrinets, and Petri nets (!) arise as particular cases of WSN in which every function is afne and isdened by matrices satisfying simple conditions: see Fig. 3. This point of view is new and putsthe decidability results of [29,7] in a clearer perspective.

    An algorithm of independent interest was developed to compute the limit of a nonnegative afnefunction (the algorithm extends to nonnegative afne functions with real coefcients).

    We leave open the questions for which ? appears in Figs. 2 and 4. It would be interesting tocomplete the global picture afforded by WSN and afne WSN by answering these questions.

    Finally, we have mentioned that well-structured transition systems capture many innitestate systems beyond generalized Petri nets (see for instance [11]). Clearly, any useful such tran-sition system ought to be dened from recursive functions over some domain, so that WSNcould in principle represent such systems as well. But there is one limitation of WSN, name-ly niteness of their dening sets of functions, which could be lifted to make WSN more ex-pressive. Are there innite state systems which would require this generality to be modelledconveniently by WSN, and if so, what would become of our Figs. 2 and 4 in this more generalcontext?

    Acknowledgments

    We thank Jrme Leroux and Philippe Schnoebelen for useful comments on earlier versions ofthis paper and we are very grateful to the anonymous referees, who read our manuscript with greatcare and who contributed numerous suggestions to clarify its presentation; in particular, we arethankful for their question concerning the -recursivity of strongly increasing functions, which wehad overlooked in an earlier proposition.

  • A. Finkel et al. / Information and Computation 195 (2004) 129 29

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